A switch is provided, which can receive a data communication at an edge of a network. The network may be made up of a plurality of switches. The switch may generate a flow channel based upon an identified source and destination for the data communication. The data communication can be routed across the plurality of switches based on minimizing a number of hops between a subset of the plurality of switches and in accordance with the flow channel.
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15. A switch fabric, comprising:
a first switch of the switch fabric to:
receiving a data packet;
determine a packet flow for the packet, wherein packets of the packet flow include same values in one or more header fields;
generate a flow channel for the packet flow, wherein the flow channel indicates a set of paths for the packet flow in the network;
select a forwarding port from a set of ports associated with the flow channel based on respective loads on the set of paths, wherein a respective port corresponds to a different path of the set of paths; and
forward the data packet via the selected forwarding port across the switch fabric; and
a second switch of the switch fabric to receive the forwarded data packet, wherein the selection of the forwarding port facilitates adaptive routing in accordance with the flow channel generated by the first switch.
1. A switch, comprising:
a processor; and
a memory unit coupled to the processor and including instructions that when executed, cause the processor to:
receive a data packet at an edge of a network, wherein the network comprises a plurality of switches;
determine a packet flow for the packet, wherein packets of the packet flow include same values in one or more header fields;
generate a flow channel for the packet flow, wherein the flow channel indicates a set of paths for the packet flow in the network;
select a forwarding port from a set of ports associated with the flow channel based on respective loads on the set of paths, wherein a respective port corresponds to a different path of the set of paths; and
forward the data packet via the selected forwarding port, wherein the selection of the forwarding port facilitates adaptive routing across a subset of the plurality of switches in accordance with the flow channel.
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16. The switch fabric of
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This application claims the benefit of U.S. Provisional Patent Application No. 62/852,273, filed May 23, 2019, entitled “Network Switch,” U.S. Provisional Patent Application No. 62/852,203, filed May 23, 2019, entitled “Network Interface Controller,” and U.S. Provisional Patent Application No. 62/852,289, filed May 23, 2019, entitled “Network Computer System,” the disclosures of which are incorporated herein by reference in their entirety for all purposes.
The invention(s) described herein were made with U.S. Government support under one or more of the contracts set forth below. The U.S. Government has certain rights in these inventions.
Contract Title
Customer/Agency
Contract Reference
FastForward-2
Lawrence Livermore
Subcontract B609229
National Security,
under prime contract
LLC/Dept of Energy
DE-AC52-07NA27344
BeePresent
Maryland
H98230-15-D-0020;
Procurement Office
Delivery Order 003
SeaBiscuit
Maryland
II98230-14-C-0758
Procurement Office
PathForward
Lawrence Livermore
Subcontract B620872
National Security,
under prime contract
LLC/Dept of Energy
DE-AC52-07NA27344
DesignForward
The Regents of the
Subcontract 7078453
University of California/
under prime contract
Dept of Energy
DE-AC02-05CII11231
DesignForward-2
The Regents of the
Subcontract 7216357
University of California/
under prime contract
Dept of Energy
DE-AC02-05CII11231
As network-enabled devices and applications become progressively more ubiquitous, various types of traffic as well as the ever-increasing network load continue to demand more performance from the underlying network architecture. For example, applications such as high-performance computing (HPC), media streaming, and Internet of Things (IOT) can generate different types of traffic with distinctive characteristics. As a result, in addition to conventional network performance metrics such as bandwidth and delay, network architects continue to face challenges such as scalability, versatility, and efficiency.
The present disclosure, in accordance with one or more various embodiments, is described in detail with reference to the following figures. The figures are provided for purposes of illustration only and merely depict typical or example embodiments.
The figures are not exhaustive and do not limit the present disclosure to the precise form disclosed.
The present disclosure describes systems and methods that can accommodate exascale computing, e.g., perform data-intensive tasks such as simulations, data analytics, artificial intelligence workloads at exascale speeds. In particular, an HPC network or interconnect fabric is provided that may be Ethernet-compatible, able to connect to third-party data storage, and can be built using a switch component that is extremely high bandwidth, e.g., on the order to 12.8 Tb/s/dir per switch with, e.g., 64 200 Gbps ports that support large network creation with very low diameter (e.g., only three network hops). Moreover, low latency can be achieved through novel congestion control mechanisms, adaptive routing, and the use of traffic classes that allows for flexibility in terms of bandwidth shaping, priority, and routing policy.
Using adaptive routing techniques, a revised routing methodology is incorporated into a network where minimal routing is further classified into preferred minimal routes and normal minimal routes across the entire path. Where the preferred route is one that will prefer a more direct minimal route, resulting in less hops to the destination.
Typically, traffic can be injected into switch fabric 100 via an ingress port of an edge switch, and leave switch fabric 100 via an egress port of another (or the same) edge switch. An ingress edge switch can group injected data packets into flows, which can be identified by flow ID's. The concept of a flow is not limited to a particular protocol or layer (such as layer-2 or layer-3 in the Open System Interface (OSI) reference model). For example, a flow can be mapped to traffic with a particular source Ethernet address, traffic between a source IP address and destination IP address, traffic corresponding to a TCP or UDP port/IP 5-tuple (source and destination IP addresses, source and destination TCP or UDP port numbers, and IP protocol number), or traffic produced by a process or thread running on an end host. In other words, a flow can be configured to map to data between any physical or logic entities. The configuration of this mapping can be done remotely or locally at the ingress edge switch.
Upon receiving injected data packets, the ingress edge switch can assign a flow ID to the flow. This flow ID can be included in a special header, which the ingress edge switch can use to encapsulate the injected packets. Furthermore, the ingress edge switch can also inspect the original header fields of an injected packet to determine the appropriate egress edge switch's address, and include this address as a destination address in the encapsulation header. Note that the flow ID can be a locally significant value specific to a link, and this value can be unique only to a particular input port on a switch. When the packet is forwarded to the next-hop switch, the packet enters another link, and the flow-ID can be updated accordingly. As the packets of a flow traverses multiple links and switches, the flow IDs corresponding to this flow can form a unique chain. That is, at every switch, before a packet leaves the switch, the packet's flow ID can be updated to a flow ID used by the outgoing link. This up-stream-to-down-stream one-to-one mapping between flow ID's can begin at the ingress edge switch and end at the egress edge switch. Because the flow ID's only need to be unique within an incoming link, a switch can accommodate a large number of flows. For example, if a flow ID is 11 bits long, an input port can support up to 2048 flows. Furthermore, the match pattern (one or more header fields of a packet) used to map to a flow can include a greater number of bits. For instance, a 32-bit long match pattern, which can include multiple fields in a packet header, can map up 2{circumflex over ( )}32 different header field patterns. If a fabric has N ingress edge ports, a total number of N*2{circumflex over ( )}32 identifiable flows can be supported.
A switch can assign every flow a separate, dedicated input queue. This configuration allows the switch to monitor and manage the level of congestion of individual flows, and prevent head-of-queue blocking which could occur if shared buffer were used for multiple flows. When a packet is delivered to the destination egress switch, the egress switch can generate and send back an acknowledgement (ACK) in the upstream direction along the same data path to the ingress edge switch. As this ACK packet traverses the same data path, the switches along the path can obtain the state information associated with the delivery of the corresponding flow by monitoring the amount of outstanding, unacknowledged data. This state information can then be used to perform flow-specific traffic management to ensure the health of the entire network and fair treatment of the flows. As explained in more detail below, this per-flow queuing, combined with flow-specific delivery acknowledgements, can allow the switch fabric to implement effective, fast, and accurate congestion control. In turn, the switch fabric can deliver traffic with significantly improved network utilization without suffering from congestion.
Flows can be set up and released dynamically, or “on the fly,” based on demand. Specifically, a flow can be set up (e.g., the flow-ID to packet header mapping is established) by an ingress edge switch when a data packet arrives at the switch and no flow ID has been previously assigned to this packet. As this packet travels through the network, flow IDs can be assigned along every switch the packet traverses, and a chain of flow IDs can be established from ingress to egress. Subsequent packets belonging to the same flow can use the same flow IDs along the data path. When packets are delivered to the destination egress switch and ACK packets are received by the switches along the data path, each switch can update its state information with respect to the amount of outstanding, unacknowledged data for this flow. When a switch's input queue for this flow is empty and there is no more unacknowledged data, the switch can release the flow ID (i.e., release this flow channel) and re-use the flow-ID for other flows. This data-driven dynamic flow setup and teardown mechanism can obviate the need for centralized flow management, and allows the network to respond quickly to traffic pattern changes.
Note that the network architecture described herein is different from software-defined networks (SDN's), which typically uses the OpenFlow protocol. In SDN, switches are configured by a central network controller, and packets are forwarded based one or more fields in the layer-2 (data link layer, such as Ethernet), layer-3 (network layer, such as IP), or layer-4 (transport layer, such as TCP or UDP) headers. In SDN such header-field lookup is performed at every switch in the network, and there is no fast flow ID-based forwarding as is done in the networks described herein. Furthermore, because the OpenFlow header-field lookup is done using ternary content-addressable memory (TCAM), the cost of such lookups can be high. Also, because the header-field mapping configuration is done by the central controller, the setup and tear-down of each mapping relationship is relatively slow and could require a fair amount of control traffic. As a result, an SDN network's response to various network situations, such as congestion, can be slow. In contrast, in the network described herein, the flows can be set up and torn down dynamically based on traffic demand; and packets can be forwarded by a fixed-length flow ID. In other words, flow channels can be data driven and managed (i.e., set up, monitored, and torn down) in a distributed manner, without the intervention of a central controller. Furthermore, the flow ID-based forwarding can reduce the amount of TCAM space used and as a result a much greater number of flows can be accommodated.
Referring to the example shown in
Note that the operations described above can be performed substantially at line speed with little buffering and delay when the first packet is received. After the first packet is processed and scheduled for transmission, subsequent packets from the same flow can be processed by switch 110 even faster because the same flow ID is used. In addition, the design of the flow channels can be such that the allocation, matching, and deallocation of flow channels can have substantially the same cost. For example, a conditional allocation of a flow channel based on a lookup match and a separate, independent deallocation of another flow channel can be performed concurrently in nearly every clock cycle. This means that generating and controlling the flow channels can add nearly no additional overhead to the regular forwarding of packets. The congestion control mechanism, on the other hand, can improve the performance of some applications by more than three orders of magnitude.
At each switch along the data path (which includes switches 110, 106, and 104), a dedicated input buffer can be provided for this flow, and the amount of transmitted but unacknowledged data can be tracked. When the first packet reaches switch 104, switch 104 can determine that the destination fabric address in the packet's fabric header matches its own address. In response, switch 104 can decapsulate the packet from the fabric header, and forward the decapsulated packet to host 116. Furthermore, switch 104 can generate an ACK packet and send this ACK packet back to switch 110. As this ACK packet traverses the same data path, switches 106 and 110 can each update their own state information for the unacknowledged data for this flow.
In general, congestion within a network can cause the network buffers to fill. When a network buffer is full, the traffic trying to pass through the buffer ideally should be slowed down or stopped. Otherwise, the buffer could overflow and packets could be dropped. In conventional networks, congestion control is typically done end-to-end at the edge. The core of the network is assumed to function only as “dumb pipes,” the main purpose of which is to forward traffic. Such network design often suffers from slow responses to congestions, because congestion information often cannot be sent to the edge devices quickly, and the resulting action taken by the edge devices cannot always be effective in removing the congestion. This slow response in turn limits the utilization of the network, because to keep the network free of congestion the network operator often needs to limit the total amount of traffic injected into the network. Furthermore, end-to-end congestion control usually is only effective provided that the network is not already congested. Once the network is heavily congested, end-to-end congestion control would not work, because the congestion notification messages can be congested themselves (unless a separate control-plane network that is different from the data-plane network is used for sending congestion control messages).
In contrast, the flow channels can prevent such congestion from growing within the switch fabric. The flow channel mechanism can recognize when a flow is experiencing some degree of congestion, and in response can slow down or stop new packets of the same flow from entering the fabric. In turn, these new packets can be buffered in a flow channel queue on the edge port and are only allowed into the fabric when packets for the same flow leave the fabric at the destination edge port. This process can limit the total buffering requirements of this flow within the fabric to an amount that would not cause the fabric buffers to become too full.
With flow channels, the switches have a reasonably accurate state information on the amount of outstanding in-transit data within the fabric. This state information can be aggregated for all the flows on an ingress edge port. This means that the total amount of data injected by an ingress edge port can be known. Consequently, the flow channel mechanism can set a limit on the total amount of data in the fabric. When all edge ports apply this limit action, the total amount of packet data in the entire fabric can be well controlled, which in turn can prevent the entire fabric from being saturated. The flow channels can also slow the progress of an individual congested flow within the fabric without slowing down other flows. This feature can keep packets away from a congestion hot spot while preventing buffers from becoming full and ensuring free buffer space for unrelated traffic.
In general, flow channels can define a path for each communication session across the switch fabric. The path and amount of data belonging to each flow can be described in a set of dynamically connecting flow tables associated with each link of the switch fabric. On every ingress port, edge and fabric, a set of flow channel queues can be defined. There can be one queue for each flow channel. As packets arrive, they either can be assigned to a flow channel on an edge port, or have been assigned to a flow channel by the link partner's egress fabric port on a fabric ingress port. The flow channel information can be used to direct the packets into the appropriate flow channel queue.
Crossbar switch 210 includes crossbars which can be configured to forward data packets and control packets (such as ACK packets) among the communication ports. EFCT logic block 212 can process packets received from an edge link and map the received packets to respective flows based on one or more header fields in the packets. In addition, EFCT logic block 212 can assemble FGFC Ethernet frames, which can be communicated to an end host to control the amount of data injected by individual processes or threads. IFCT logic block 214 can include the IFCT, and perform various flow control methods in response to control packets, such as endpoint-congestion-notification ACKs and fabric-link credit-based flow control ACKs. OFCT logic block 216 can include a memory unit that stores the OFCT and communicate with another switch's IFCT logic block to update a packet's flow ID when the packet is forwarded to a next-hop switch.
In one embodiment, switch 202 is an application-specific integrated circuit (ASIC) that can provide 64 network ports that can operate at either 100 Gbps or 200 Gbps for an aggregate throughput of 12.8 Tbps. Each network edge port may be able to support IEEE 802.3 Ethernet, and Optimized-IP based protocols as well as Portals, an enhanced frame format that provides support for higher rates of small messages. Ethernet frames can be bridged based on their L2 address or they can be routed based on their L3 (1Pv4//1Pv6) address. Optimized-IP frames may only have an L3 (1Pv4/1Pv6) header, and are routed. Specialized NIC support can be used for the Portals enhanced frame format, and can map directly onto the fabric format of network 100, e.g., a fabric format that provides certain control and status fields to support a multi-chip fabric when switches/switch chips, such as switches 102, 104, 106, 108, and 110 are connected and communicate with each other. As alluded to above, a congestion control mechanism based on flow channels can be used by such switches, and can also achieve high transmission rates for small packets (e.g., more than 1.2 billion packets per second per port) to accommodate the needs of HPC applications.
Switch 202 can provide system-wide Quality of Service (QoS) classes, along with the ability to control how network bandwidth is allocated to different classes of traffic, and to different classes of applications, where a single privileged application may access more than one class of traffic. Where there is contention for network bandwidth, arbiters select packets to forward based on their traffic class and the credits available to that class. Network 100 can support minimum and maximum bandwidths for each traffic class. If a class does not use its minimum bandwidth, other classes may use the unused bandwidth, but no class can get more than its maximum allocated bandwidth. The ability to manage bandwidth provides the opportunity to dedicate network resources, as well as CPUs and memory bandwidth to a particular application.
In addition to support for QoS classes, switch 202 effectuates flow channel-based congestion control, and can reduce the number of network hops, e.g., in a network having a dragonfly topology, from five network hops to three. The design of switch 202, described in greater detail below, can reduce network cost and power consumption, and may further facilitate use of innovative adaptive routing algorithms that improve application performance. A fabric created by a plurality of switches, such as a plurality of switches 202 may also be used in constructing Fat-Tree networks, for example when building a storage subsystem for integration with third-party networks and software. Further still, the use of switch 202 enables fine-grain adaptive routing while maintaining ordered packet delivery. In some embodiments, switch 202 may be configured to send the header of a packet from an input port to an output port before the full data payload arrives, thereby allowing output port load metrics to reflect future loads, thereby improving adaptive routing decisions made by switch 202.
By providing flow-specific input queues, the switch can allow each flow to move independently of all other flows. The switch can avoid the head-of-queue blocking problem, which is common with shared input buffers. The flow-specific input queue also allows the packets within a single flow to be kept in order. When a flow passes through the switches, a flow-specific input queue on each input port can be allocated for this flow and these input queues become linked, effectively forming one long queue that reaches across the entire fabric for this flow, and the packets of this flow can be kept in order.
The progress of successful delivery of packets belonging to a flow can be reported by a sequence of ACKs generated by the edge port of an egress switch. The ACK packets can travel in the reverse direction along the data path traversed by the data packets and can be forwarded by the switches according to the forwarding information maintained in flow tables. As ACK packets travel upstream, they can be processed by each switch's input queue manager, which can update the corresponding flow's state information based on information carried by the ACK packets. The ACK packets can have a type field to provide advanced information about the downstream data path, such as congestions. A switch's input queue manager can use this information to make decisions, such as throttling the transmission rate or changing the forwarding path, about the pending data packets currently buffered in its input queues. In addition, the input queue manager can update the information carried in an ACK packet based on state information of a buffered flow, so that the upstream switches can make proper decisions. For example, if an input queue for a given flow is experiencing congestion (e.g., the amount of data in the queue is above a predetermined threshold), the input queue manager can update an ACK packet that is being forwarded to the next upstream switch to include this congestion information.
If an ACK corresponds to the last packet of a flow, a switch can determine that there is no more unacknowledged data for that flow. Correspondingly, the switch can free the flow channel by removing the corresponding entry in the flow table.
As mentioned above, the input queue manager at each switch can maintain information about transmitted but unacknowledged data of a given flow.
In the example in
Correspondingly, because the input queue for the given flow at switch 262 has six queued data units, and two additional data units are in transit between switches 252 and 262, the flow_extent value at switch 252 is 9. Similarly, the flow_extent value at switch 202 is 13, because there are three data units stored in the input queue at switch 252 and one data unit in transit between switches 202 and 252.
In general, a flow channel can remain allocated to a single flow until all the ACKs for all the packets sent on the flow channel have been returned. This means that flow channel table entries can remain active for longer near the fabric ingress edge port than near the egress edge port. If a single packet is injected into the network, a flow channel can be allocated for the ingress edge port and then another flow channel can be allocated for the next fabric link the packet traverses and so on, until the last flow channel is allocated when the packet reaches the last fabric link. Each allocation can generate a flow ID, denoted as variable “flow_id,” to identify the entries of the flow tables of the fabric link. (More details on flow channel tables are provided in the description below in conjunction with
At the input queue of each switch, the flow channel table entries can indicate each flow's state information, including the flow_extent value, from this point downstream to the flow's egress destination edge port. Packets received on the local input port can increase this flow_extent value by the amount of incoming data, and ACKs can reduce the flow_extent by the amount of acknowledged, delivered data.
When a packet reaches the final destination egress port, an ACK packet can be generated and returned for that packet. This ACK can be routed using the data path information stored in the corresponding entry of the flow channel tables at every switch along the data path. Optionally, the ACK packet itself does not need to carry path information and therefore can be small and light weight. If no other data packet is sent on the flow, the ACK can release each flow channel in the reverse order. Once released, the flow channel at each switch can be allocated to a different flow.
If another packet follows the first packet on the same flow, the ACK corresponding to the second packet would need to be received before the flow channel can be released at a given switch. In one embodiment, the flow channel can only be released when ACKs for all the transmitted packets of the same flow have been returned.
Typically, various protocols may require in-order packet delivery. The flow channels can be used to guarantee this delivery order, even when the fabric uses adaptive routing for load balancing across multiple data paths. If packets between an ingress edge port and an egress edge port, perhaps in a different switch on the far side of the fabric, are injected at a very low rate, then each packet injected could reach its destination and return an ACK back to the source before the next packet is injected. In this case, each packet can be a lead packet and free to take any path across the fabric, using the best available dynamic adaptive routing choice. This is possible because the first packet can define the flow's path through the fabric.
Now assume that the packet injection rate is increased slightly to the point where the next packet of the same flow is injected before the current packet's ACK has returned to the source. The second packet can pass the ACK of the first packet somewhere along the flow's data path. Beyond this passing point, the ACK will have released the flow channels allocated to the first packet, because the flow_extent value associated with the first packet is returned to zero when the ACK is processed by the flow channel's logic. Meanwhile, the second packet can now define a new flow, because it is again causing flow channels to be allocated on each of the subsequent fabric links. This second packet, while it is causing flow channels to be allocated beyond the passing point, can be forwarded to a different path based on dynamic adaptive routing. On the other hand, before the passing point, the second packet can extend the outstanding flow created by the first packet to include the second packet. This means the first packet's ACK may not reduce the flow_extent value to zero and the flow channels may remain active before the passing point. It also means that the second packet may follow the exact path taken by the first packet up to the passing point. Note that while it is following the previous packet, the second packet cannot arrive at the egress edge port before the first packet does, and therefore correct packet order can be maintained.
If the injection rate for this flow is increased further, the second packet will pass the first packet's ACK at a location closer to the destination edge port. It is also possible that a third, fourth, fifth, or additional packet may enter the fabric before the first packet's ACK is returned to the source edge port, depending on the data packet injection rate of this flow and the data packet-ACK round trip delay. The maximum packet rate can depend on the size of the packets and the bandwidth of the links. The round trip delay of the data packet and ACK can be an important parameter for a fabric implementation and can be used along with the maximum packet rate to calculate the maximum required number of flow channels for each link. Ideally, a design can provide a reasonable number of unallocated flow channels regardless of the traffic pattern. The demand for the number of flow channel can be high when a large number of packets arriving at an ingress edge port have different destinations and these packets have small sizes and high injection rates. In the most extreme case, each packet could be allocated a different flow channel. These flow channels are freed when the packets' ACKs are returned. Correspondingly, the number of flow channels needed can be calculated as ((Packet rate)*(Average packet to ACK round trip latency)).
Note that packet rate on a single flow channel is not to be confused with packet rate on a link. If the traffic pattern is such that many small packets are being sent to different destinations, then successive packets sent onto the link can have different destinations. This means that each packet could belong to a different flow and could be the only packet to use the corresponding flow channel. In this example, the link can experience a high packet rate, but the packet rate of individual flows can be low. Optionally, a number of ACKs (e.g., 48 ACKs) can be aggregated together into a single ACK frame for transmission over a link and protected by a Frame Check Sequence (e.g., a 32-bit FCS). For example, the ACKs can occupy 25 bits each, and there can be a 9-byte overhead to the frame. That is, the overhead per ACK on a full size frame is approximately 9/(25/8*48)*100%=6%. The logic can optimize the number of ACKs per frame so an ACK does not need to wait too long to be aggregated when the ACKs are arriving slowly. For example, the ACK aggregation logic block can use three timers to manage ACK transmission based on the activity of an outgoing link. These timers can be started when a new ACK arrives at the ACK aggregation logic block. If the outgoing link is idle, a first timer, which can for example be set at 30 ns, can be used to hold the ACK while waiting for additional ACKs to arrive. When this timer expires, all the ACK received within the corresponding time window can be aggregated into one frame and transmitted onto the outgoing link. If the outgoing link is busy, a second timer, which can for example be set at 60 ns, can be used to wait for additional ACKs. Using this second timer can allow more ACKs to be aggregated into a single frame, and this frame can be transmitted only if a predetermined number of ACKs are collected. Note that due to the Ethernet framing constrains, some numbers of ACKs in a single frame can use less wire bandwidth per ACKs than other numbers of ACKs. If no efficient number of ACKs are collected, and the outgoing link remains busy sending normal data packets, then a third timer, which can for example be set at 90 ns, can be used. Once this third timer expires, all the ACKs that have been collected can be aggregated in a frame and transmitted onto the link. By using these three timers, the system can significantly reduce the overhead of sending ACKs on the outgoing link.
In some examples, the ingress edge port of a switch can encapsulate a received data packet with a fabric header, which allows the packet to be forwarded using flow channels.
Crossbar switch 210 may comprise separate, distributed crossbars routing data/data elements between input and output ports. In some embodiments, and as illustrated in
Request crossbar 210a is used to send requests from an input to a targeted output age queue. Grant crossbar 210b is used to return a grant back to the input to satisfy a request. In particular, grant crossbar 210b returns a pointer indicating where a packet is within an input buffer. It should be noted that a grant is returned when there is space in the output for the corresponding packet. Grant crossbar 201b may also optionally return a credit for requested space in the output. It should be noted that grants are returned when there is a landing spot for a packet at the output, e.g., an output port 220c, so packets cannot be blocked (though they can face transient contention for resources).
It should be understood that in accordance with various embodiments, a credit protocol may be used to guarantee that there is a landing space for a request at the output. Accordingly, a credit crossbar 210c may be used to return credit for requested space in the output.
A data crossbar 210d is used to move granted packets from an input buffer to a targeted output buffer. An Ack crossbar 210e is used to propagate Ack packets from output ports 220c to input ports 220b. Acks are steered in accordance with a state kept in an output flow channel table.
It should be understood that data crossbar 210d moves multi-clock packets with both headers and data, while the other four crossbars (request crossbar 210a, grant crossbar 210b, credit crossbar 210c, and Ack crossbar 210e) move only single-clock packet headers. All five crossbars use the same architecture with row buses and column buses within an 8×4 matrix of 32 dual-port tiles (as described below).
Referring back to
Each crossbar switch 210 has sixteen inputs 220b, one for each port in its row, and eight outputs 220c, one for each port in its column. Row buses can be driven from each source in a row to all eight crossbars in that row (one-to-all). Arbitration can be performed at the crossbar from the sixteen row buses in that row to the eight column buses in a given column. Buffering can be provided at each 16×8 crossbar for each of the row buses in order to absorb packets during times when there is contention for a column bus. In some embodiments, a non-jumbo packet is kept off a row bus unless there is room for the entire packet in the targeted crossbar input buffer. Due to area constraints, jumbo packets are allowed to go even if there is not sufficient space (crossbar input buffer only sized to sink a non-jumbo packet) with the row bus being blocked until the packet wins arbitration and space is freed as it is moved onto a column bus.
Column buses are driven from a given crossbar to each destination port within a column (all-to-all). Each destination may have another level of arbitration between the column buses from the four rows. With sixteen row buses driving eight crossbars, each feeding eight column buses, there is a 4× speedup between rows and columns. Each row has identical connections with the one-to-all row bus connections for a single row shown in row buses. Each tile will have a one (request, grant, credit) or a two (data, ack) clock delay per tile depending on the crossbar. This gives a maximum seven or fourteen clock delay to get between the leftmost and rightmost columns. Credit returns routed through credit crossbar 210c may have a one clock delay per tile, and therefore, can take a maximum of seven clocks to complete transmission.
It should be noted that each column may have identical connections with the all-to-all column bus connections for a single column, and there may be a two clock delay per tile, resulting in a six clock delay to get from the top row to the bottom row. It should also be understood that both row and column buses both use the aforementioned credit-based protocol to determine when they are able to send. In the case of row buses, the source port maintains credit counts for the input buffers of the crossbars within that row. For the data crossbar, care is needed to determine when a packet is allowed to go on a row bus. If grants targeting a particular crossbar input buffer all go through a single queue, space for the packet at the head of the queue is required before starting the packet transfer. If the grants are distributed across multiple queues, in order to prevent small packets from locking out large packets, a packet transfer does not start unless there is space for an entire max sized packet in the buffer. In this way, once a packet transfer on a row bus starts, it will not stop until the entire packet has been transferred. Accordingly, crossbar input buffers are configured to be large enough to handle the maximum packet size plus additional space to cover the worst case round trip (packet send to credit return). This will not be the case for jumbo packets. To save on buffering area, the crossbar input buffers are only deep enough to handle a non-jumbo sized MTU (1500 bytes) with a jumbo packet being allowed to block a row bus while waiting to gain access to the targeted column bus.
For column buses, each crossbar maintains credit counts for the input buffers at each destination port in that column. Unlike row buses, there is no requirement that credits be available for a maximum-sized packet before starting transfer of that packet on a column bus. Individual words of the packet will move as credits become available. Therefore, the input buffer at the destination for each column bus needs to only be big enough to cover the worst case round trip (packet to credit).
A fair round-robin arbitration may be used between the sixteen sources for each destination. For the data crossbar 210d, once a source wins arbitration, it keeps control of the destination column bus until the entire packet has been sent. Each output grants a limited amount of packet payload so it is expected that contention for a given column bus should be fairly limited when larger packets are involved. Because of this, a round-robin arbitration is expected to be sufficient even with possibly large differences in packet size among requesters.
Parts of switch 202 associated with output functions generally operate on frames within the switch fabric format, and have a fabric header, even, for example, for a frame arriving and leaning on an Ethernet port within a single switch 202.
Age queue output control [INVENTOR QUESTION—WHERE IS “
Requests come into the output control block via a column bus from each row of matrix 30. Each column bus feeds an independent FIFO (e.g., first-in-first-out shift register or buffer) with space in the FIFO managed via credits. The FIFOs may be sized (24 deep) to cover a round-trip plus additional space to allow requests to be moved out of the crossbars 210a-210e and prevent head-of-line blocking. Prior to writing into a FIFO, a request may be checked for a valid error correcting code (ECC). If the ECC check has either a multi bit error (MBE) or a single bit error (SBE) in the destination field (i.e. it has been routed to the wrong port), the request is considered to be an invalid request, and is discarded with an error being flagged.
Least recently used (LRU) arbitration may be performed between column bus FIFOs to choose which FIFO gets forwarded to age queue management. As requests are removed from each FIFO, credits are returned to the corresponding crossbar. The row with which an incoming column bus corresponds can be dependent both on where in the matrix the tile is located, and as well as which half of the tile the block is in.
The output buffer (OBUF) makes requests to the output control block for sending reduction and maintenance packets across a link. These requests may be given the highest priority. A FIFO with 8 locations can be used to buffer these reduction/maintenance packet requests while they wait for resources. Reduction packets need not use flow channels, and maintenance packets may use loopback to create a flow so that checking for flow channel availability or flowing through the OFCT to create a grant is not needed. Reduction and maintenance packets also need not use any space in the output buffer so that no check of space is required. Rather, a check for the link partner input butter may be performed. If allowed, a shaping queue (SQ) or virtual channel (VC) can be granted, blocking any grants from the age queue path from being granted during that cycle.
The size of the next request to be processed from the output buffer is checked against max_frame_size. If it exceeds this setting, the request is not processed and an error flag is set. This will result in the output buffer request path being blocked until a warm reset is performed. The error flag will stay set until the reset is done. The condition can also be released by increasing the setting of max_frame_size to a value above the size of the stuck output buffer request. The size used in the comparison may be the size indicated in the output buffer request (which may include a 4-byte frame checksum (FCS) used on the wire).
Each input may be given the same fixed allocation of age queue space. This age queue space is large enough to reserve a location for each SQ/VC with enough additional space to cover a request/credit round-trip. It is up to the input to manage the space it is given across its SQs/VCs. This allocation (fixed_al/oc) is programmable via a control and status register (CSR) in each input queue (INQ), and can be, e.g., in the range of 64-96 locations. The remaining age queue space (8K−64*fixed_al/oc) may be shared space that is available to all inputs. The shared space can be managed by the output with it moving incoming requests from static to shared space as they arrive if there is room in the shared space, subject to per-input limits. When moving a request to the shared space, a credit is returned, e.g., immediately, via credit crossbar 210c, with the request marked in the age queue as being in the shared space.
When a request is granted, if it is marked as using the shared space, the shared space is credited. If it is not marked as using shared space, the request is considered to have used the static space, and a credit is returned to the input with the grant.
Due to conflicts in credit crossbar 210c, it is possible that credits may not be sent every clock period. Accordingly, a FIFO provides buffering for these transient disruptions. Space in this FIFO is required before taking a request from the request crossbar. A FIFO with a depth of 32 locations can be used to limit the chances of it ever backing up into request crossbar 210a. The shared space may have limits for how much space any input (from an input port 220b) can take. These limits can be set as a percentage of the available space. For instance, if the limit is set to 50%, if one input port is active, it has access to 50% of the buffer space, with two active input ports, each gets 37.5% ((space_used_by_I+pace_left*0.5)/2=(50%+50%*0.5)/2), with three active input ports, each gets 29.2% ((space_used_by_2+space_left*0.5)/3=(75%+25%*0.5)/3), and so on. Additionally, the total space used by the active input ports can be limited to the given total (50%, 75%, 87.5%). Thus, the space allocated to each of input port 220b may vary dynamically by how many inputs ports are currently active. The addition of an active input port causes other active inputs ports to give up their space which is then taken by the new input.
Given that division is not something easily done in hardware, the aforementioned age queue credit management function can be implemented as a lookup table 310 with 64 entries 312. The number of inputs currently active in the age queues 320 indexes 315 the lookup table 310. The values 314 in the lookup table 310 reflect the limit of the number of shared space locations any input can take along with the total space they can consume as a whole. Thus, it is up to software to program the values 314 in the lookup table 310 according to how much total shared space there is and what percentage each input port is allowed to take. As more input ports 220b become active, each input port 220b is allowed less space, and the total space available increases. Incoming requests from input ports 220b that are above this limit, or in total, exceed the total space limit, are not allowed to take more shared space. In order to track the number of active input ports 220b in the age queues, a set of 64 counters 316 (one for each input port) is used. These count up when a request is put in the age queues 320 and count down as they are taken out (i.e., granted). A count of the number of non-zero counts 319 is used as an index into the lookup table 310. In addition, in order to manage the shared space, an additional set of 64 counters 318 may be used to track the current usage of the shared space by each input. There may also be a single counter 334 that can be used to track overall shared space usage. These counters are compared against the current quotas to determine if a request is allowed to use the shared space or not. Counters 316, 318 can be, e.g., 13 bits wide, to provide sufficient coverage of the maximum value of an object that may be somewhat less than 8K.
Age queues 320 may use a single storage RAM 321 that has 8K locations in it. These locations can be dynamically allocated to 32 separate queues (one for each SQ/VC) with each consisting of a linked-list of locations within the storage RAM 321. This gives each SQ/VC the ability to take more space as needed.
An age queue 320 can be created with a front pointer 322 pointing to the front of the queue, and a next pointer 324 for each location pointing the next item in the queue. The last location in the queue may be indicated by a back pointer 326. Items are taken from the front of the queue and inserted at the back of the queue. In addition to the above data structures, each queue has a FIFO 328 of entries at its head. These FIFOs 328 may ensure that a queue can sustain a request every clock with a multi-clock read access time from the request RAM 321. When a new request arrives, if the head FIFO 328 328 for that queue is not full, it bypasses the request RAM 321, and can be written directly into the head FIFO 328. Once requests for a given age queue are being written to the request RAM 321, subsequent requests are also written to the request RAM 321 to maintain order. The bypass path can be used again once there are no more requests for that age queue in the request RAM 321, and there is room are the head FIFO 328. When a request is read from a head FIFO 328, and there are corresponding requests queued in the request RAM 321, a dequeue is initiated. One head FIFO 328 may be read at a time, such that a single dequeue operation can be initiated each clock period. Logic may be included to handle the various race conditions between an ongoing or imminent enqueue operation and a head FIFO 328 being read.
The aforementioned ECC protection used in the age queue RAM 321 can be extended to the FIFOs 328 to protect the data path flops. The resulting structure may include 8K flops (32 queues×5 deep×SQ-bits wide). When generating the ECC, the age queue number can be included in the calculation (but not stored) as an extra check of the free list management. When the ECC is checked, the request can be considered to be in error if there is an MBE or there is an SBE in the queue number bits.
A free list RAM can be a simple FIFO which is initialized with pointers to all 8K entries whenever a reset is performed. A count can be maintained to keep track of how may entries are valid within the free list. As entries are taken, they are popped off the front of the FIFO and used. As entries are returned, they are pushed onto the back of the FIFO. Some number of entries, e.g., three entries,) at the head of the free list can be kept in flops so they are available for quick access. As with the head FIFOs for the age queues, ECC is carried through the flops to provide protection. The resulting structure may have minimal flops (57=3 deep×19-bits wide).
In order to support full performance for small packets, age queues support both an enqueue and a dequeue every clock period. The operations across the data structures for an enqueue operation are discussed below, and can differ depending on whether the queue being written is empty or not.
In some cases, a simultaneous enqueue and dequeue to a specific queue is easily handled as they are using and updating separate fields. Some specialized scenarios may arise, e.g., when a dequeue operation empties the age queue. In order to handle this scenario, a dequeue occurs first logically, followed by an enqueue operation. Accordingly, an empty flag is seen as being set when the queue is emptied by the dequeue operation, and then cleared due to the enqueue operation.
The arbitration alluded to above can be performed among requests that are permitted to be granted subject to input buffer management, output buffer management, and flow channel quotas. Arbitration can also be halted if there are no credits for the OFCT input FIFO. In some embodiments, arbitration may be performed at two levels. First, traffic shaping arbitration can be used to arbitrate between the SQs. A Deficit Round-robin arbitration can be used to arbitrate between VCs within a given SQ. Traffic shaping arbitration may use a series of token buckets to control the bandwidth of each SQ as follows: eight leaf buckets, one for each SQ; four branch buckets; and a single head bucket.
Arbitration can be divided into three groups with a first group having the highest priority, followed by a second group, which in turn is followed by a third group. For the first and second groups, arbitration may be handled in the same way among eligible SQs. A ×8 round-robin arbitration can be performed between the SQs for each of the eight priority levels (eight parallel round-robin arbitrations). A fixed arbitration can be performed between priority levels. For example, group 3 arbitration has no priorities, and therefore is simply a single ×8 round-robin arbitration.
For arbitration in the first group, the priority for each comes from the setting in the leaf buckets. For arbitration in the second group, priority comes from the setting in the branches of the leaf buckets. In all cases, the buckets which are checked to be eligible for that group, are also the buckets from which packet size tokens are obtained if that request wins arbitration.
Regarding age queue 320 selection, packets can be classified in order to select the SQ to which their request is forwarded. This allows traffic associated with an application to be shaped differently from traffic originating from a different application or a different traffic class. This can be useful on the edge ports which connect to a NIC in that the applications will have been configured to use a share of the resources on the node, and similarly will be granted a proportion of the network bandwidth. In accordance with one embodiment, this classification is performed by classifying the packets into a traffic class identifier (FTAG), e.g., a 4-bit code that is part of the fabric frame header, and a VLAN ID (VNI) as the packet ingresses into the fabric. The FTAG and VNI may then be used as the packet egresses the fabric to select the shaping queue.
A configuration register can be used to map FTAGs to SQs. This configuration matches the corresponding configuration in the in queue. When the output buffer requests or returns link partner credits, it converts a given FTAG to an SQ. For packet injection, the FTAG is found in R_TF_OBUF_CFG_PFG_TX_CTRL. For test generation, the FTAG is found in the test control register. When the reduction engine (RED) requests a credit return, the FTAG is found in ret_cdtltag. When a reduction frame is removed from the output stream and link partner credits need to be returned, the FTAG is found in the frame header.
Regarding the SQs discussed herein, each age queue 320 may have 32 SQs that are addressed by {SQ, VC}. The 3-bit SQ 330 can be considered a shaping function, and the VC selects one of four queues within that shaping function. For Ethernet egress (edge) ports, the VC is not needed for deadlock avoidance. Accordingly, all 32 SQs 330 can be available. In such a scenario, the SQ 330 can be selected by adding the SQ base from R_TF_OBUF_CFG_FTAG_SQ_MAP to the lower bits of the VNI. The 5-bit sum defines the {SQ,VC} to send to the age queue. It should be noted that when injecting frames on an egress port, a VNI is not available, and therefore, an SQ base can be directly used. For fabric links, the SQ 330 is taken from the upper three bits of the SQ base. The VC can be taken from the frame header when returning credits for reduction frames, or from the appropriate control CSR (R_TF_OBUF_CFG_TEST_CTRL or R_TF_OBUF_CFG_PFG_TX_CTRL) when injecting frames.
A link partner input buffer management can depend on the type of device to which the link is attached. Devices such as switch 202 may use credit-based flow control where each credit represents a cell of storage in the input buffer. Other devices may use standard Ethernet pause or priority pause-based flow control. Requests which are marked to terminate locally (lac term set) need not consider link partner input buffer flow control and need not update any associated counters. Link partner space need not be considered when the link is in the draining state.
For credit-based flow control, the link partner input buffer can be divided into eight buffer classes. Each SQ 330 can be assigned to one of these 8 buffer classes. Credits are maintained for each of the buffer classes with each credit representing 32 bytes of storage in the link partner input buffer. In order to allow credit-based flow control to work with various devices (switch, enhanced NIC), each of which may have different cell sizes, the cell size is a programmable value in units of 32 bytes.
There may be two sets of VCs with each SQ 330 assigned to one set. A maximum frame size worth of space can be reserved for each VC, and each VC set can have a different maximum frame size. The remainder of the link partner input buffer is shared dynamic space usable by any SQ/VC, subject to per VC and buffer class limits.
The size that comes with the request represents the size of the packet on the wire which includes a 4-byte FCS. This gets converted to an internal 2-byte FCS at the link partner before writing the packet to the link partner input buffer so the crediting needs to account for this difference, which can be a factor at the boundary of the cell size. For instance, for a 96 byte cell, a size that is 97 or 98 will take a single cell. In order to know when this happens, the request includes a correction term which is calculated as: req.len_correct=(byte_len % 16)==1 or 2.
Further validation of this term is required to convert it to whatever the cell size boundary may be. It will be valid when the length just exceeds the cell size. With this, the validated fen_correct term can be determined by: len_correct=(((16−byte size) % (2*32−byte cell size))==1) & req. len correct
An example of how these values work for a few cell and packet sizes is illustrated in the table below:
Length Correct Calculation
Size
Req
Size
Cell Size
len_correct
Credit
(bytes)
len_correct
(16 B units)
(32 B units)
Modulo result
len_correct
Taken
64
0
4
2
0
0
2
65
1
5
2
1
1
2
66
1
5
2
1
1
2
67
0
5
2
1
0
3
96
0
6
3
0
0
3
97
1
7
3
1
1
3
9B
1
7
3
1
1
3
99
0
7
3
1
0
4
128
0
8
4
0
0
4
129
1
9
4
1
1
4
130
1
9
4
1
1
4
131
0
9
4
1
0
5
The size that comes with the request uses 8-byte units and the link partner input buffer cell size is a multiple of 32 bytes (32*y where y=cell size from CSR). First, the 8−byte size is converted to a 16−byte size (ROUNDUP((8−byte size)/2)). Also, the cell size is converted to 16 byte units (2*y). Mathematically, the number of cells a request will use can be calculated by: ROUNDDN(((16−byte size)+2*y−1−len_correct)/(2*y))=# of cells
While a divide operation is possible in hardware, due to timing reasons, a divide operation cannot be done in the critical path of the arbitration. Instead, an alternate credit management is used. That is, credits are maintained in units of 32 bytes. When a request wins arbitration, the number of credits taken is adjusted by the maximum error term (2*y−1) using the calculation: ROUNDDN(((16-byte size)+2*y−1)/2)=Maximum 32 byte credits needed. Because this calculation overestimates the credit required for the packet, on the following clock, a modulo operation (X=(16-byte size) MOD 2*y, y=32-byte cell size from CSR) can be performed to determine what the actual remainder is. This value along with the len_correct term are used to adjust the credit counter. The formula used to create the adjustment value (adf_val) for Xis: If (X==0) adj_val=y−1 else if (X==1 and fen_correct) adj_val=y else adj_val=ROUNDDN((X−1)/2)
The table below illustrates a request credit example for 96 byte cells showing the values used across several packet lengths for the 96 byte cells of the switch input buffer (y=3).
Request Credit Example for 96-Byte Cells
Packet Size
Packet Size
Credit
Modulo
Corrected
(bytes)
(16-byte units)
Taken
Result
len_correct
adj_val
Credit Taken
48
3
4
3
0
1
3
64
4
4
4
0
1
3
80
5
5
5
0
2
3
96
6
5
0
0
2
3
97
7
6
1
1
3
3
98
7
6
1
1
3
3
99
7
6
1
0
0
6
128
8
6
2
0
0
6
If a request is filtered before being forwarded to the link partner input buffer, the output buffer logic returns the SQ and VC so they can be used to return the credits to the appropriate credit counters. No size is required since the packet size is always the same, the length of a reduction frame (69-byte or 16-byte size=5).
The local (master) side of the link maintains a count of the number of packets sent from each VC across both sets (8 total), a count of amount of packet (in 3 2-byte quantities) sent to each VC (4), and a count of the amount of packet (in 32-byte quantities) sent for each buffer class(8). The link partner (slave) side of the link maintains the same set of counts with them being sent over the link periodically. The difference between the master and slave counts is a count of the number of packets in the link partner input buffer from each VC across both sets and a count of the amount of space (in 32-byte quantities) currently occupied by each VC and each buffer class. A count is also maintained of the total amount of space used across all packets. A summary of the counters is as follows: master_vcx_cnt[4]/slave_vcx_cnt[4]—master and slave counts of the number of packets sent to each VC in set X; master_vcy_cnt[4]/slave_vcy_cnt[4]—master and slave counts of the number of packets sent to each VC in set Y; master_bc_cnt[8]/slave_bc_cnt[8]—master and slave counts of the amount of space occupied by each buffer class in units of 32-bytes; master_vc_cnt[4]/slave_vc_cnt[4]—master and slave counts of the amount of space occupied by each VC in units of32-bytes; master-tot-cnt/slave-tot-cnt—master and slave counts of the total amount of space occupied in units of32-bytes.
All counters are set to zero on a warm reset. They are also forced to zero when the link is in the draining state or when the DBG_RESET CSR bit to clear their state is set. The output buffer filter will steer a reduction packet to something other than the path to the link partner input buffer. In this case, a signal can be returned along with the SQ and VC of the packet. Again, the length is not required as the size of these packets is fixed. This information is used to adjust the appropriate master credit counts.
A request is allowed to participate in arbitration if either its VC count is 0 (indicating its one statically assigned slot is available) or there is space for a max sized frame in the dynamic space (subject to the targeted buffer class and VC limits). There can be a single programmable value for max frame size which is used across all VCs and SQs. The request validation for input buffer space can be addressed using credit-based flow control.
Credit-based flow control can be used to divide a dynamic space in two ways, each independent of each other: first, based on a limit of how much dynamic space each of the four VCs can take; and second, based on a limit to how much dynamic space each of the eight buffer classes can take. In both cases, the limits are set as a percentage of the available space. For a given packet, space should be made available in both its targeted VC and buffer class. For instance, if each space has its limit set to 50%, if one is active, it has access to 50% of the buffer space, with two active, each space gets 37.5% ((50+50*0.5)/2), with three active, each space gets 29.2% ((75+25*0.5)/3), and so on. Also, the total space used by those spaces that are active can be limited to the given total (50%, 75%, 87.5%). Accordingly, the space allocated to each varies dynamically by how many are currently active. When an additional one goes active it causes others that are active to give up some of their space which is then taken by the new one.
Like the division function discussed above, this function is implemented as a lookup table. For the VC space in this example, there are 16 entries with each entry specifying the space available to each VC along with the total space available across all VCs. For the buffer classes, there may be 256 entries with each entry specifying the space available to each buffer class along with the total space available across all buffer classes. Space for each is expressed in 2048-byte units. The depth of each table is sufficient to cover all combinations of active members (VCs or buffer classes), with each being able to have an independent setting for their percentages. With this, it is up to software to program the values in the table according to how much total dynamic space there is and what percentage each is allowed to take across all possible combinations. As more become active, each is allowed less space and the total available increases. Requests for spaces that are above this limit, or in total above the total limit, are not allowed to take more dynamic space.
A VC or buffer class is considered active either if it has a request in an age queue, or if it has outstanding credits for link partner input buffer space. As an example, consider there are only 4 spaces (16 entry table) with percentages set as SPACE0(50%), SPACE1(40%), SPACE2(30%), SPACE3(10%), and a total dynamic space of 16 KB. This results in the values, in quantities of 16-bytes presented in the buffer space example table below.
Buffer Space Example
Index
SPACE3
SPACE2
SPACE1
SPACE0
Total
0
N/A
N/A
N/A
N/A
N/A
1
N/A
N/A
N/A
512
512
2
N/A
N/A
410
N/A
410
3
N/A
N/A
319
398
717
4
N/A
307
N/A
N/A
307
5
N/A
250
N/A
416
666
6
N/A
255
339
N/A
594
7
N/A
202
270
337
809
8
102
N/A
N/A
N/A
102
9
94
N/A
N/A
469
563
10
94
N/A
377
N/A
471
11
75
N/A
299
374
748
12
95
284
N/A
N/A
379
13
78
234
N/A
389
701
14
80
239
319
N/A
638
15
79
236
315
394
1024
As an example, the values in the row for index 7 are calculated as: Total %=0.5+(1−0.5)*0.4+(1−0.5−(1−0.5)*0.4)*0.3=0.79; SPACEO=(0.5/(0.5+0.4+0.3))*0.79*1024=337; SPACEI=(0.4/(0.5+0.4+0.3))*0.79*1024=270; SPACE2=(0.3/(0.5+0.4+0.3))*0.79*1024=202; Total=337+270+202=809
As noted above, and referring back to
Routing in the switch fabric may be controlled by a fabric routing function (FRF) implemented in switch 202. An example FRF component 400 is illustrated in
In particular, FRF component 400 determines a preferred port with preferred ports discriminator 402B to forward each frame presented at the input port based on: a received frame's destination fabric address (DFA); the frame's current routing state (where the frame is along its path, and the path(s) it took to reach its current routing state); the switch fabric routing algorithm and configuration; and load metrics associated with the output port (the aforementioned preferred port to which the frame is to be forwarded) using busy ports filters.
FRF component 400 may include a routing algorithm table 408 that may be embodied as a software configurable table that determines valid choices based on the frame's current routing state. Valid choices are decisions such as whether a local minimal, global minimal, local non-minimal, or global non-minimal path is allowed to be chosen for the frame's next hop. The routing state includes information such as the VC the frame was received on, and whether it is in the source, the destination, or an intermediate group. The routing algorithm table 408, along with the adaptive selection function or logic 412A (described below), also determines the VC to be used for the frame's next hop.
Frame routing with unicast DFAs will be described as an example. However, it should be noted that the DFA of the routing request can either be in unicast of multicast format. The unicast format can include a 9-bit global ID field (global_id), a 5-bit switch ID field (switch_id), and a 6-bit endpoint ID field (endpoint_id). The global ID can uniquely identify a group within the network. Specifically, it identifies the final group to which the frame must be delivered. The switch ID uniquely identifies a switch within the group identified by global ID. The endpoint ID field, together with the global ID and switch ID identify the endpoint, connected to the edge of the network fabric, to which the frame is to be delivered. This field is mapped to a port or set of ports on the switch identified by global ID and switch ID.
The multicast format includes a 13-bit multicast ID field (multicast_id). This field is mapped by FRF component 400 to a set of ports on the current switch to which the frame is to be forwarded.
From this information, FRF component 400 determines an updated routing state for the frame, which is then carried within the frame. For example, to effectuate routing in a dragonfly topology, a frame's current state may be gleaned from the frame's VC (discussed above). Based on algorithmic switch fabric routing rules specified for the switch fabric (the selection of which is described below), FRF component 400 determines a particular VC to be used for the frame's next hop to avoid any deadlocks. Additional routing state information can be provided depending on where the frame is along its path, e.g., whether the frame is in its source group, in an intermediate group, or in its destination group. It should be noted that FRF component 400 performs port filtering (described in greater detail below) using permitted ports filter, operational ports filter, busy ports filters, etc. to determine if a preferred port to which a frame is to be forwarded is currently faulty, busy, absent, etc.
Switch 202 distributes load information between switches. The FRF component 400 receives the load measurement of and from its associated output port. The FRF component 400 receives summary load information from its associated input port for a neighboring switch. Each FRF component 400 exchanges load information with all other FRF instances within the same switch. FRF component 400 provides summary load information to its associated output port for delivery to a neighboring switch. Through the load distribution mechanism, each FRF component 400 learns the load measured at each output port of its switch. As well, each FRF learns the summary load information for all neighboring switches.
It should be noted that FRF component 400 can support frame multicasting. When a multicast DFA is received, FRF component 400 determines a set of ports to which the frame associated with the multicast DFA should be forwarded. The set of ports can be determined by accessing a lookup table that maps software-configured multicast fabric addresses to output ports. This helps avoid problems associated with duplicate multicast frame copies.
Filtering may be applied to the three categories of candidate ports, e.g., operational port filtering, useable port filtering, and busy port filtering. Port filtering as applied herein can be used to reduce the set of valid ports considered as path candidate ports by identifying and removing absent and/or faulty ports from consideration.
Operational port filtering (or non-operational port filtering) can refer to the removal of non-operational ports from sets of ports being considered as candidates for routing, e.g., preferred minimal path candidate ports 502, non-preferred minimal path candidate ports 504, and non-minimal path candidate ports 506. That is, switch 202 may identify certain ports as being non-operational. These non-operational ports may be reported in a non-operational port mask. It should be noted that in some embodiments, software can force additional ports of switch 202 to be considered as non-operational using a non-operational port CSR, e.g., when a port(s) is to become disconnected as a result of planned maintenance.
Usable (or unusable) port filtering may involve filtering out candidate ports for consideration that would normally have been acceptable, but due to faults within network 100, for example, the candidate ports have become unacceptable/unusable for reaching one or more destination switches, destination groups (of switches), etc., but remain acceptable or usable for reaching one or more other destination switches. In some embodiments, global fault table 406A can be used to block global minimal path port candidates and global non-minimal path port candidates depending on a destination group of the frame being routed. For example, candidate ports that lead to an intermediate group (of switches) without connectivity to a particular destination group (of switches) can be excluded from consideration when routing frames to that destination group, although the same candidate ports may not necessarily be blocked out for other destination groups. The global fault table 406A can be determined or indexed by the global_id field of the frame's DFA.
In some embodiments, an exception list maintained by exception list table 404A may be used to conditionally exclude port candidates depending on the destination group or switch to which the frame is being routed. It should be noted that that exception list table 440A may be used to identify preferred global minimum path ports. Accordingly, use of the exception list table 404A to exclude port candidates is done when it is not being used to identify preferred global minimum path ports.
It should be noted that knowledge regarding which ports are busy in a neighboring switch can be used to determine if the ports that connect to a neighboring switch are poor candidates to receive a forwarded frame based on whether the neighboring switch will subsequently need to forward the frame to a port that is already busy. For example, when considering candidate ports for use in global minimal routing, the ports connected to a neighboring switch are poor candidates if the neighboring switch's global ports that connect to the frame's destination group are all busy. Similarly, when in the destination group and considering candidate ports for use in local nonminimal routing, the ports connected to a neighboring switch are poor candidates if the neighboring switch's local ports that connect to the frame's destination switch are all busy.
Accordingly, busy port filtering can be performed by FRF component 400 by using busy port masks to remove heavily loaded ports from being considered candidate ports. It should be noted that in some embodiments, heavily loaded ports are removed from consideration when other candidate ports that are not heavily loaded exist. Otherwise, when non-heavily loaded ports do not exist, busy port filtering will not remove the heavily loaded ports from consideration. FRF component 400 maintains four masks of busy ports, that is, ports whose load exceeds a software-defined threshold: local switch busy port mask; global non-minimal busy global port mask; global non-minimal busy local port mask; remote switch busy port mask. Information from these masks is communicated between switches to populate the remote switch.
A local switch busy port mask can be applied to minimal path candidate ports as well as to local non-minimal path candidate ports. The FRF generates a 64-bit Is_busy_port_mask by comparing each port's local_load to a software defined threshold. Ports with loads higher than this threshold are marked as busy in this mask.
A global non-minimal busy port global mask can be applied to global ports of global non-minimal path candidate ports. The FRF generates a 64-bit gnmbgp_mask by comparing each port's gnmgp_load to a software defined threshold. Ports with loads higher than this threshold are marked as busy in this mask.
A global non-minimal busy local port mask can be applied to local ports of global non-minimal path candidate ports. The FRF generates a 64-bit gnmblp_mask by comparing each port's gnmlp_load to a software defined threshold. Ports with loads higher than this threshold are marked as busy in this mask.
A destination group dependent busy-port mask, obtained from a remote switch busy global port table can be applied to global minimal path candidate ports. Correspondingly, when a frame is being routed in its destination group, a destination switch-dependent busy-port mask, obtained from a remote switch busy local port table can be applied to local non-minimal path candidate ports.
Upon applying the aforementioned filtering or down-selection stage, a set of surviving path candidate ports 508 can result. That is, a reduced number of candidate ports can be determined after removing non-operational and unusable port candidates, heavily loaded, a set of path candidate ports remains. In some embodiments, a pseudo-random selection process is used to further reduce the number of surviving path candidate ports 508 to a determined threshold number of ports associated with each category of candidate ports (preferred minimal path candidate ports, non-preferred minimal path candidate ports, and non-minimal path candidate ports). In some embodiments, that threshold number of candidate ports may be four candidate ports per category. If the minimum threshold number of candidate ports is not met, no candidate ports from that category are removed from consideration.
In some embodiments, this pseudo-random selection (or down selection) of candidate ports can be weighted. That is, weights, e.g., weights between 0 and 15, can be assigned to each port per the CSR configuration. This weighting can be used to influence the probability with which individual candidate ports are chosen such that higher-weighted ports have a great chance of being chosen. For example, a weight of 15 results in a port having 15 times greater likelihood of being selected in the pseudo-random selection process. In some embodiments, candidate ports may be filtered into four groups (GW1, GW2, GW4, GW8) based on their assigned weights, where a candidate port can belong to multiple groups depending on the assigned weight (e.g., a candidate port with weight 1 belongs only a one weight group, while a candidate port with weight 5 belongs to two groups (GW1 and GW4, i.e., 1+4=5), and a candidate port with weight 15 belongs to all four groups (Gw1, GW2, GW4, GW8, i.e., 1+2+4+8=15). The number of candidate ports in each group can be determined (nW1, nW2, nW4, nW8), and pseudo-random selection is applied to each group to select one candidate port from each group (cW1, cW2, cW4, cW8). The weight of each group can be computed, along with their total weight: wW1=nW1; wW2=2*nW2; wW4=4*nW4; wW8=8*nW8; wtotal=wW1+wW2+wW4+wW8. A fifth pseudo-random selection can be performed to choose a number j in the range O . . . Wtotal−1. One of the candidates, cW1, cW2, cW4, cW8 is chosen as the down selected candidate port based on the value of j as follows: If j<wW1, choose cW 1; else If j<wW1+wW2, choose cW2; else If j<wW1+wW2+wW4, choose cW4; else choose cW8.
The number of ports in each group can vary from request to request due to changing operational status and load, and may also be dependent on the configuration of the global fault table 406A. It may also vary depending on the type of port at which the routing request is being performed, i.e., edge port versus local port, for example. On the assumption that operational status and load do not change too quickly, and that the configuration of global fault table 406A should not vary greatly for different global_id values, each pseudo-random generator's previously generated value is simply used as the offset for biasing its next value. If an offset value is out of range (>n−1), it is brought within range through truncation of upper bits. mgnm=4 candidate ports are produced by the pseudo-random down selection process. Each candidate can be produced through a separate copy of the weighted pseudo-random down selection logic 410C, described above.
It should be noted that the same candidate port may be chosen by more than one of the instances/iterations of the weighted pseudo-random down selection logic 410C, in effect reducing the number of candidates ports that are chosen. The probability of the same candidate port being chosen by more than one of the mgnm=4 global non-minimal, weighted pseudo-random selectors decreases with increasing numbers of candidates ports to choose from. In the context of a dragonfly topology, for example, and a network with full global bandwidth, at an edge port in the source group there are potentially about 48 possible global non-minimal candidates ports: 16 global ports and 32 local ports. If a local hop has been taken, the next hop is a global hop, thereby reducing the number of candidate ports to about 16. However, if the network is tapered such that it only supports one quarter of full global bandwidth, there may only be 4 global candidates to choose from after a local hop has been taken. The probability of selecting four unique global non-minimal candidate ports shows the probability of four unique candidates ports being chosen for varying numbers of possible candidates to choose from. The probability of selecting n unique global non-minimal candidate ports shows the probability of varying numbers of the four selected candidates being unique when there are only 16, 8, or 4 candidate ports to choose from.
FRF component 400 can use received remote switch busy port masks to generate the aforementioned Remote Switch Busy Global Port Table of busy port masks that identify ports connected to neighboring switches that should be avoided as global minimal path candidates based on the destination group to which the frame is being routed. Similarly, the received remote switch busy port masks can also be used to generate the aforementioned remote switch busy local port table of busy port masks that identify ports connected to neighboring switches that should be avoided as local non-minimal path candidates, when routing in the destination group, based on the destination switch to which the frame is being routed.
The rs_busy_port_masks are used in assessing the suitability of neighboring switches for reaching specific destination groups via global minimal paths and specific destination switches via local nonminimal paths or via local minimal paths. Each FRF instance corresponding to a local port or a global port can be configured to generate a 64-bit rs_busy_port_mask. The generated mask is delivered to the partner switch connected to that port. Similarly, the partner switch can also generate and return an rs_busy_port_mask.
Each FRF instance communicates the rs_busy_port_mask, that it received from its partner switch, to all other FRF instances in the switch using the port status ring (which connects the tiles of a switch and communicates status and load information amongst the ports on the switch). Each FRF instance captures all rs_busy_port_masks such that all FRF instances learn the remote busy port status being provided by all neighboring switches. Each FRF instance uses the rs_busy_port_masks that it receives to generate the busy port tables described in the Remote Switch Busy Global Port (RSBGP) Table and Remote Switch Busy Local Port (RSBLP) Table.
Generation of the rs_busy_port_mask is a two-step process. The first step is to compare each port's local_load to a software configurable generating an intermediate mask of all ports that are individually busy. This intermediate mask is formed as the status of each port is received from the port status ring interface. Ports that are classified as non-operational are also recorded as busy in the intermediate mask. The second step takes link bundling into account such that a port is only marked as busy in the rs_busy_port_mask if it and all other ports, that are part of the same bundle, are marked as busy in the intermediate mask. Either all ports that are members of the same bundle are marked as busy in the rs_busy_port_mask, or none are. Global ports that are part of the same bundle all connect to the same remote group. Local ports that are part of the same bundle all connect to the same remote switch within the current group.
As the rs_busy_port_masks are used to determine whether the switch that generated the mask is a good candidate for routing a frame to another group or to another switch in the current group, bundling is used to provide a consistent view of the generating switch's suitability when the busy status across its links, that connect to the destination group or to the destination switch in the current group, is inconsistent. The rationale for the treatment of bundling described here is that the switch generating the rs_busy_port_mask remains a candidate for reaching the destination group, or the destination switch in the current group as long as it has at least one link to the destination group or switch that is not busy; adaptive routing at the switch that generated the rs_busy port_mask should direct the frame to the non-busy link.
Ports must be included in either the bundled ports mask CSR or the unbundled ports mask CSR (both of which are part of the static description of the wiring) in order for them to be marked as busy in the rs_busy port_mask. The second step is performed as each frame is received from the port status ring. The bundled port masks are scanned to identify the bundles and the ports that they contain. In addition, the unbundled port mask is consulted to identify any other ports, that are not members of a bundle, but whose busy status should also be included in the generated rs_busy_port_mask.
A different software-defined threshold is used in computing the rs_busy_port_mask because of the larger latency involved in communicating and processing the rs_busy_port_mask and in delivering a frame subject to this mask to the remote switch that generated the mask. Because of the larger latency, it may be useful to require a port to be more loaded before it is considered to be so busy that it is not a good candidate for receiving additional frames from a remote switch. A busy remote port should be sufficiently loaded such that it remains loaded throughout the time it would take to receive a frame that has been subject to the mask.
The aforementioned RSBGP table stores busy port masks indexed by destination group (global_id). Again, the RSBGP table is used in the evaluation of global minimal paths consisting of a hop to a neighboring switch which has a global port connected to the destination group to filter out ports of the current switch which are poor choices for use in reaching the destination group because the corresponding global port or ports, of the neighboring switch reached by the filtered-out ports of the current switch, are too heavily loaded.
The RSBLP table stores busy port masks indexed by destination switch (switch_id), and again, can be used in the evaluation of local non-minimal paths consisting of a local hop to a neighboring switch followed by another local hop to the destination switch. For topologies, such as fat-tree, where a local minimal path can consist of a local hop to a neighboring switch followed by another local hop to the destination switch, the RSBLP Table can also be used in the evaluation of local minimal paths. The RSBLP table is used to filter out ports on the current switch which are poor choices for use in indirectly reaching the destination switch because the neighboring switch's port or ports that connect to the destination switch are too heavily loaded.
It should be noted that the RSBGP table and the RSBLP table are never both accessed for the same routing request. The former is accessed when routing a frame that is not in the destination group, and the latter is accessed only when routing a frame that is in the destination group. Therefore, both are implemented within the same memory, termed the Remote Switch Busy Port
Provided that there is at least one valid candidate port, the various busy port filters (Busy Ports Filter, Local Non-Minimal (LN) Busy Port Filter, Global Non-Minimal (GN) Busy Port Filter) may not be allowed to collectively block all candidate ports. If there are viable candidate port choices, they are allowed, despite being “poor” choices, if there are no better choices. Otherwise, an empty route response will be generated for the routing request when routes are actually available.
To prevent an incorrect empty route response from being generated, the first stage of the preferred and non-preferred minimal path busy port filters (
It should be noted that there will be no minimal path candidate ports if minimal routing is disabled (Permitted Ports Filter). There will be local non-minimal path candidate ports only if local non-minimal routing is enabled (Candidate Local Non-Minimal Path Ports). There will be global non-minimal path candidates only if global non-minimal routing is enabled (Candidate Global Non-Minimal Path Ports). Local and global non-minimal routing are generally not both enabled simultaneously. When the first stage of the preferred and non-preferred minimal path busy port filters and the local non-minimal busy port filter are disabled due to the conditions described above, the only candidate ports that will be seen at an adaptive selection stage (described below) will be poor candidates because they will all be ports that lead to other switches whose ports (that connect to the destination group or to the destination switch) are heavily loaded. This is because these are the only candidate ports that were being blocked by the filters that are being disabled and, without these filters being disabled, there are no other candidates.
The adaptive selection stage will choose between these remaining/surviving candidate ports, which are all poor, based on their biased local loads (Local Load and Load Value Selection), although their local loads will not necessarily reflect the reason why they are poor. Their poor character can be the result of high downstream load on certain ports of the other switches reached by these candidate ports. It is because the adaptive selection stage may not be able to see how poor these candidates are that the coordination between the different busy port filters, described herein may be used. If each busy port filter decides independently whether or not to disable its RSBGP Table and RSBLP Table-based filters, situations such as the following could occur. The non-preferred minimal path busy port filter might produce one or more candidates, which are not poor, without any of its filter stages disabled. The preferred minimal path busy port filter might only be able to produce one or more candidate ports by disabling both of its filter stages. Thus, all of the candidate ports it is able to produce are poor. At the adaptive select stage, the down-selected, not poor, non-preferred minimal path candidate ports, are compared against the down-selected, poor, preferred minimal path candidate ports. However, the adaptive selection stage lacks visibility into how poor the preferred minimal path candidates are, so it may select a poor preferred minimal path candidate over a not poor non-preferred minimal path candidate.
An alternative to the busy port filter coordination mechanism described herein, would be for all of the busy port filters to act independently, but for the minimal path and the local non-minimal path busy port filters to each forward a signal through to the adaptive select stage to indicate if their respective candidate ports are poor choices due to busy ports at downstream switches. If they are, the adaptive selection function may de-prioritize their candidate ports in favor of other ports. The result would be the same as is achieved by the coordination, described herein, between the various busy port filters.
As illustrated in
In some embodiments, preferred minimal path candidate ports are preferentially selected over non-preferred minimal path candidate ports, and minimal path candidate ports are to be preferentially selected over non-minimal path candidate ports. To accomplish this preferential selection, a bias value can be added to each candidate port's load before the adaptive selection comparison is performed. The bias value used can be configured using CSRs, and can vary depending on the type of path to which it is being applied (i.e., non-preferred minimal, preferred minimal, an non-minimal), the traffic class of the frame being routed, and where the frame is along its path. For example, frames belonging to a low-latency traffic class can be more strongly biased towards minimal paths versus frames in other traffic classes to have a greater likelihood of achieving/comporting with low-latency requirements or needs. Frames may also be increasingly biased towards minimal path routes the close the frames are to their destination.
In particular, load values represent the busyness of switch 202's ports and are used in evaluating the load-based port masks and in comparing candidate ports during the adaptive route selection process. Load-based port masks are used in the busy port filters to remove ports that are poor candidates, based on current load, from the set of candidates being considered. There are a number of different types of load values used within Switch and some are communicated to neighboring Switch devices. These load values are described below, and illustrated in
Several load metrics are computed and used in determining which port to route a frame when there is more than one port to which the frame can be routed. The load metrics are also used in generating busy-port masks, which as described above, are used to remove heavily loaded ports from consideration.
There are five load metrics described herein: local load; group load; global non-minimal global port load; mean global load; and global non-minimal local port load.
Regarding local load, the load of each of a switch's output ports (e.g., output ports 220c of switch 202) is continuously being evaluated and provided to the corresponding FRF instance as the 8-bit value local_load. Larger values represent higher load. The current load present at each output port is measured by the output control age queue block. The output port load is provided by each age queue instance to the FRF instance (of FRF component 400) that is associated with the input side of the same port. The load value provided to the FRF is an 8-bit value termed the local_load. The age queue determines the local load based on a combination of the amount of traffic enqueued, waiting to go out that port and the amount of traffic enqueued at the opposite side of the link in the link-partner Switch device's input buffer. The calculation and configuration of local_load is performed later. Each FRF instance distributes the local_load value it receives from its associated age queue instance to all other FRF instances. In this way, each FRF instance learns the current local_load of every output port.
When the port loads of candidate ports are being compared to determine the best port to route a frame to, it is the port's local_load value that is used for ports being considered for minimal path routing, and for ports being considered for local non-minimal path routing.
Group load is a measure of how suitable a dragonfly group is for use as the intermediate group in a global non-minimal path. The 8-bit group load value is not computed by a switch, such as switch 202, but is software-configurable. Software might use a measure of the network injection load present across the input side of the group's edge ports in deriving the group_load value, or might determine the group_load value based on a policy of discouraging use of certain groups as non-minimal intermediate groups, perhaps based on the jobs or services that are running in the groups. That is, the group_load value is intended to be representative of the amount of local traffic within a group.
A network management stack sets the group_load value by periodically writing to a CSR. The software-configured group load value is communicated across global links. FRF instances associated with global links forward the group_load value that they receive from their link-partner in the group at the opposite side of the link, to all other FRF instances in the switch. In this way, each FRF instance learns the group load values of the groups at the opposite end of each of the global links terminated by the switch.
In terms of global non-minimal global port load (gnmgp_load), this load is a metric used in assessing a global port's suitability for use in directing a frame to the intermediate group reached by the global link connected to the global port. The gnmgp_load is nominally equal to the maximum of the global port's local_load and the group_load value being received from the group reached by the global link. However, through field G-NMGP_EN_GRP LD in CSRR_TF_FRF_CFG-_LOAD_CTRL, the group_load component can be excluded.
When the port loads of candidate ports are being compared to determine the best choice port to route a frame to, it is the port's gnmgp_load value that is used for global ports being considered for global non-minimal path routing.
Mean global load (mean_global_load) is intended for use in assessing a switch's suitability for use in reaching any intermediate group that is directly connected to that switch. The mean_global_load value is an 8-bit value equal to the arithmetic mean of the gnmgp_load values of all of the switch's global ports. Ports that are classified as non-operational by either hardware or software are excluded from the calculation. The ports to include in the mean_global_load computation are determined from CSR R TF FRF CFG GNM global ports.
For any port whose load is included in the mean_global_load computation, if a group_load value is not being received for that port from the port status ring, either because link partner data is not being received for that port or because the link partner data that is being received is not global link data, the contribution of that port to the mean_global_load is based solely on that port's local load. It should be understood that the aforementioned port status ring communicates status and load information amongst the ports on a switch (e.g., input and output ports 220b and 220c, respectively, of switch 202). The computed mean_global_load value is communicated across local links. FRF instances associated with local links forward the mean_global_load value, that they receive from their link-partner in the local switch at the opposite side of the link, to all other FRF instances in a switch. In this way, each FRF instance learns the mean_global_load values of the local switches at the opposite end of each of the local links terminated by the switch, and can use these values in global non-minimal path selection.
Global non-minimal local port load metrics are computed by each FRF instance, distributed between ports in a switch using the port status ring, and distributed between switches. The global non-minimal local port load is a metric for use in assessing a local port's suitability for use in directing a frame to an intermediate group of a global non-minimal path. The global non-minimal local port load takes into account the load on the local port as well as the suitability of the local group switch, to which the port connects, for use in reaching an intermediate group. A port's gnmlp_load value is equal to the maximum of the port's local_load and the mean_global_load reported by the port's partner switch. Through software configuration it is possible to remove the mean_global_load component such that a port's gnmlp_load becomes simply equal to its local_load.
The gnmlp_load value is an 8-bit value. The computed gnmlp_load value is based on the local_load and mean_global_load values distributed via the port status ring. Each FRF instance computes the gnmlp_load value for all of its switch's ports at which link partner data for a local link is being received. If load status information is not being received from a port's partner switch, the gnmlp_load value for that port is set equal to the port's local_load.
For ports at which link partner data for a global link is being received, the port's gnmlp_load value is set to the value computed for gnmgp_load. This is a side-effect of an implementation optimization in which the same storage is used for gnmgp_load and gnmlp_load since, for any given port, at most one of the two is valid. The global non-minimal local port load metric is not used on ports where global link partner data is being received.
The fabric routing process described in the preceding subsections is performed for every frame received. Switch determines whether to perform packet-by-packet adaptive routing (using this value) or flow based adaptive routing (where this value is used for the first packet in each flow) according to the ordering requirements of the traffic.
As has been described herein, a switch, such as switch 202 supports minimal and non-minimal path routing. It should be understood that minimal paths are based on the destination. If a destination NIC is local, an output port that connects to the destination switch is selected. If the destination is in another group, the packet is routed to a switch within the local group that is connected to the destination group. In a large system, there may only be one such path, but in a small system there are likely to be many, some connected to the input switch and others to switches elsewhere within its group. The input switch selects between them.
As has been described herein, a switch, such as switch 202 supports minimal and non-minimal path routing in a network. In some embodiments, as alluded to above, the network may have a dragonfly topology, Dragonfly routing is hierarchical, distinguishing between local destinations (those in same group as source) and global destinations. Thus in a dragonfly network, a switch routes to a destination group and then to a switch within that group using two tables rather than to individual destinations using one large table.
It should be understood that minimal paths are based on the destination. If a destination NIC is local, an output port that connects to the destination switch is selected. If the destination is in another group, the packet is routed to a switch within the local group that is connected to the destination group. In a large system, there may only be one such path, but in a small system there are likely to be many, some connected to the input switch and others to switches elsewhere within its group. The input switch selects between them.
The candidate ports considered for minimal path routing are further divided into preferred and nonpreferred sets of ports where ports in the preferred set may lead to a path containing fewer hops. Non-minimal paths route packets via an intermediary switch, referred to as a root switch. Root switches are selected on a packet-by-packet basis or a flow-by-flow basis depending on the ordering requirements of the traffic.
The candidate ports considered for minimal path routing are further divided into preferred and nonpreferred sets of ports where ports in the preferred set may lead to a path containing fewer hops. Non-minimal paths route packets via an intermediary switch, referred to as a root switch. Root switches are selected on a packet-by-packet basis or a flow-by-flow basis depending on the ordering requirements of the traffic.
Non-minimal traffic is routed “up” to the root switch, and then minimally “down” to the destination. In some embodiments, intermediate root switches are selected at random so as to distribute load uniformly. The network, e.g., network 100, provides control over intermediate group selection enabling traffic to be routed towards intermediate groups that are known to be lightly loaded or away from those that have specific function or are known to be heavily loaded. Root switches may be distributed over all groups, where a non-minimal path may detect a root switch in the source group, the destination group, or any intermediate group. Global non-minimal routes take an indirect path through a root switch in an intermediate group. These paths require two global hops, one from the source group to an intermediate group, and one from the intermediate group to the destination group. Global non-minimal paths require up to three local hops, one in each group. The maximum path length is five switch-to-switch hops, whatever the system size.
Minimal routing is to be preferred as the paths are shorter, and hence the load on the network is lower. However, minimal routing alone will result in poor performance on some traffic patterns, for example when all nodes in one group communicate with nodes in one other group. Achieving good performance across a wide range of traffic patterns requires a mix of minimal and non-minimal routing.
At each hop along a frame's path, the routing modes that may be used to advance the frame along its next hop are controlled by the configuration of the FRF routing algorithm table 408. When a frame is received at a switch input port, the types of paths along with the frame may be forwarded is determined: local minimal, global minimal, local non-minimal, and global non-minimal. The set of output ports to which the frame may be forwarded is determined by the type of paths allowed at that point.
The types of paths that are allowed to be taken depends on where the frame is at along its journey between its ingress and egress ports of the network. The path types are as follows. Local minimal paths select links that are directly connected to the frame's destination switch and may be used when the frame is in its destination group. Global minimal paths may be used when the frame is not in its destination group and select either global links that directly connect to the frame's destination group or local links that connect to a switch that has working global links that directly connect to the frame's destination group. Local non-minimal paths may be used when the frame is in its destination group, or when the frame is in an intermediate group. Local non-minimal paths select local links connected to other switches in the group without regard for the frame's destination. When in the destination group, it must be possible to reach the frame's destination switch within one more hop after taking the local non-minimal hop. When in an intermediate group, it must be possible to reach a switch with a working global link that connects to the frame's destination group within one more hop after taking the local non-minimal hop. Local links, that connect to switches from which this in not possible, must not be selected.
Global non-minimal-paths may be used when the frame is in its source group and is not in its destination group. Global non-minimal-paths select either global links connected to other groups or local links connected to other switches in the source group without regard for the frame's destination. Global links must only be selected if they connect to a group that has working links connecting it to the frame's destination group. Similarly, local links must only be selected if they connect to switches that have global links that are, themselves, valid global non-minimal path choices.
Adaptive routing selects between minimal and non-minimal paths (described above) according to their current load.
In terms of minimal routing, when in the destination group, but not at the destination switch, local minimal routes are generated by looking up the switch_id field of the destination fabric address in a local switch minimal table (
Global minimal routes are generated by looking up the global_id field of the destination fabric address in a global minimal table (
Local minimal paths take at most one switch-to-switch hop, from the source switch to the destination switch, both of which are within the same group. There can be several such paths. Local non-minimal paths take two switch-to-switch hops, from the source switch to an intermediate switch, known as the Root switch, and from there to the destination switch. There are many such paths.
Global minimal paths take one global hop from the source group to the destination group. There is at most, one local hop in each of the source and destination groups. Global minimal paths require a maximum of three switch-to-switch hops whatever the system size.
In certain system configuration in which there are multiple global links connecting a source group to a destination group, a bias can arise such that the proportion of traffic injected at the source group that is distributed to each of the global links is not equal. As an example, consider the case of switch A, B, and C, all in the group X, with switch B having three global links connecting it to group V and switch C having two global links connecting it to group V. If traffic injected at switch A, destined for group V, is equally distributed between switch B and C, each of the two global links of switch C will be more heavily loaded than each of the three global links of switch B.
To enable the bias to be counteracted, the Global Minimal Table can be divided into several blocks, each of which is capable of generating a valid set of global routing choices for any destination group. On a frame by frame basis, the block that is used to service the request is pseudo-randomly selected by the FRF. Within each Global Minimal Table block instance, only a subset of the possible candidate ports that can be used to reach the destination group are populated. The subset is chosen in such a way so as to counteract the bias. The populated subsets can be varied across the different block instances such that all possible candidate ports are able to be used.
A global minimal path between an edge port in one group and an edge port in another group can require one, two, or three hops across fabric links. One hop if the switch containing the ingress edge port in the source group has a global link directly connecting it to the switch containing the egress edge port in the destination group. Two hops if the frame traverses a global link between the two groups that is connected directly to either the ingress switch in the source group or to the egress switch in the destination group. In this case, one hop across a local link in either the source group or the destination group is also required. Lastly, three hops are required if the frame traverses a global link that is not directly connected to either the ingress switch in the source group or to the egress switch in the destination group. In this case, a hop across a local link is also required in both the source and destination groups. The local link hop in the source group takes the frame from the ingress switch to the source group switch that is connected to the global link. The local link hop in the destination group takes the frame from the destination group switch that is connected to the global link to the egress switch.
When identifying minimal path candidate ports, the FRF is able to classify the candidates into a set of preferred ports and a set of non-preferred ports. The preferred ports are those that allow a global minimal path requiring two, or fewer fabric link hops. The non-preferred ports are all of the minimal path candidates that are not classified as preferred. Use of preferred paths, when available, and when not too heavily loaded reduces the average load on the system's local fabric links as it reduces the average number of local fabric links traversed per frame. Use of a preferred path may also reduce the end-to-end fabric latency experienced by the frame.
When performing local non-minimal routing, any local link can be a candidate. However, some local links may need to be excluded from consideration if they lead to a switch from which it may not be possible to reach the destination because of link or switch failures that exist within the system. A CSR controls which ports are candidates for local non-minimal path routing.
When performing global non-minimal routing, generally any global link can be a candidate. Additionally, generally any local link that reaches a switch with operational global links can also be a candidate. However, some links may need to be blocked from consideration if they lead to a switch or to a group from which it may not be possible to reach the destination group because of link or switch failures that exist within the system. CSRs control which ports are candidates for global nonminimal path routing.
When selecting a candidate port to use for global non-minimal routing, if all candidate ports, global and local, are equally likely to be selected, for many system configurations globally non-minimal traffic will not be evenly distributed among the global links leaving the group. Consider, for example, the situation of three switches, A, B, and C within a group, where each switch is connected to each other switch by four local links, and switch A and switch B each terminate 14 global links and switch C terminates 16 global links. Ingress traffic arriving at an edge port of switch A may be routed globally non-minimally to any of the global or local links terminated by switch A. If it is equally likely to be routed to any of these links, then each global link terminated by switch B will receive only 4/14th (4 local links reaching 14 global links) of the traffic routed to each global link terminated by switch A. Similarly, each global link terminated by switch C will receive 4/16th (4 local links reaching 16 global links).
To counteract this potential bias in the distribution of globally non-minimal traffic among a group's global links, when the set of global non-minimal candidate ports is being pseudo-randomly down selected to the small number of ports that will participate in the adaptive routing stage, a weighting can be applied to each of the candidate ports such that some will be more likely than others to survive the down-selection process.
Adaptive routing selects between these minimal and non-minimal paths based on their load. Adaptively routed traffic starts on a minimal path, diverting to a non-minimal path if the load on the minimal path is high (this is known as progressive adaptive routing). Such paths are said to have diverged.
A non-minimal path can be selected at the injection point or at the exit router in a source group. A local non-minimal route may be taken within the source group, an intermediate group, or a destination group if it is selected as the intermediate. Dragonfly routing algorithms allow a non-minimal path in both the intermediate and destination groups (consider a case where all traffic incoming on the global ports of a particular router is destined for NICs on another router in the group). In general however, non-minimal traffic is sufficiently well distributed as to avoid this happening, but an additional hop in the destination group may still be beneficial in cases where there is an error on a local link. Having arrived at the intermediate group the packet may take either a minimal route to the destination group or a local non-minimal route to a switch with a path to the destination group.
Again, this decision is made based on load. Having taken a hop within the intermediate group the packet must detect root and take a minimal path to the destination. Adaptive decisions are made based on load and a bias towards preferred, minimal, or non-minimal. A routing algorithm, as described above, increases the bias towards minimal paths the closer a packet is to its destination. This algorithm prefers a direct path across the intermediate group provided the load is low.
Restricted routing is used at points other than injection and root detection to prevent packets from flowing back in the direction from which they came. In a switch, cases in which a packet has taken one hop from points of injection and root detection are detected and ensured that a global port is taken. For local minimal routing, having taken a hop from the root, the packet will arrive at a switch that is connected to the destination NIC. For global minimal routing, having taken a hop from the point of injection, the packet will arrive at a switch with a global link, which must be taken. In the intermediate group, packets are allowed to take a local hop at the point of injection and having detected root. Having taken this hop, the packet will arrive at a switch that is connected to the destination group, this path must be taken.
When passing information from switch to switch, it is necessary to make instantaneous decisions about the next hop on the path. Decisions are taken using information derived from local state and information that is communicated from neighboring switches. Use of information from many different sources allows for more accurate/effective decisions. This includes information from neighbors.
Prior systems carried information on average load from switch to switch. That said, more detailed information from related or neighboring switches is more helpful. In the current switch ASIC, a set of values can include information indicative of the status of output ports of related/neighboring switches. By passing this set of values, much better routing decisions can be made. In one example, a flag is passed back from neighboring switches, with the flag having one bit for each output port. For example, with a switch having 64 outputs, a 64 bit flag would be transmitted. This is much more accurate than simply passing a global average for neighboring ports.
In accordance with one embodiment, and as described above, port busyness can be taken into account. That is, and referring back to
Filtering may be applied to the three categories of candidate ports, e.g., operational port filtering, useable port filtering, and busy port filtering. Port filtering as applied herein can be used to reduce the set of valid ports considered as path candidate ports by identifying and removing absent and/or faulty ports from consideration.
In the example of
Here, a local switch busy port mask can be applied to minimal path candidate ports as well as to local non-minimal path candidate ports, e.g., output ports of intermediate switches 704 and 706. FRF component 400 generates a 64-bit Is_busy_port_mask by comparing each port's local_load to a software defined threshold. Ports with loads higher than this threshold are marked as busy in this mask. In this example, output ports from intermediate switch 704 directed to destination switch 708 meet or surpass the software defined threshold specified for the switch fabric. Therefore, the Is_busy_port_mask indicates that the link from intermediate switch 704 to destination switch 708 is busy. On the other hand, the output ports of intermediate switch 706 to destination switch 708 are not marked or flagged with the Is_busy_port_mask as the busyness threshold is not met/surpassed, reflecting that the link between intermediate switch 706 and destination switch 708 is quiet, and not busy.
Based on the busy port mask indication that the link between intermediate switch 704 is busy, the output port of intermediate switch 704 linked to an input port of destination switch 708 is deemed to be a poor candidate, and removed from consideration. Accordingly, and in contrast to the scenario of
At operation 800, a data communication is received at an edge of a network. As described above, a switch, e.g., switch 202 may have a plurality of ports (64 ports in some embodiments), some of which may be input ports (220B), some of which may be output ports (220C). Packets may be routed between input and output ports of one switch to another one or more switches (each having their own input and output ports. Routing of packets between switches can be performed in accordance with certain characteristics of the switches that the packets may traverse along a path from a source switch to a destination switch. In some embodiments a switch 202 may be a switch at the edge of the network or switch fabric, and may receive data, e.g., data packets or frames from another switch, from a NIC, etc., either within or external to the switch fabric via an ingress port of such an edge switch.
At operation 802, a flow channel based on an identified source and destination for the data communication is generated. As described above, upon receiving injected data packets, an ingress edge switch can assign a flow ID to the flow, and the flow ID can be a locally significant value specific to a link and to a particular input port. As the packet is forward to a next hop switch, the packet enters another link and the flow ID is updated such that as the packet traverses links and/or switches, a unique chain us formed by the flow IDs corresponding to this particular flow. Subsequent packets belonging to the same flow can use the same flow IDs along the path. When packets are delivered to the destination egress switch and ACK packets are received by the switches along the data path, each switch can update its state information with respect to the amount of outstanding, unacknowledged data for this flow. When a switch's input queue for this flow is empty and there is no more unacknowledged data, the switch can release the flow ID (i.e., release this flow channel) and re-use the flow-ID for other flows. This data-driven dynamic flow setup and teardown mechanism can obviate the need for centralized flow management, and allows the network to respond quickly to traffic pattern changes.
At operation 804, the data communication is routed across a plurality of switches based on minimizing a number of hops between a subset of the plurality of switches and in accordance with the flow channel. At each hop along a frame's path, the routing modes that may be used to advance the frame along its next hop are controlled by the configuration of the FRF routing algorithm table 408. When a frame is received at a switch input port, the types of paths along with the frame may be forwarded is determined: local minimal, global minimal, local non-minimal, and global non-minimal. The set of output ports to which the frame may be forwarded is determined by the type of paths allowed at that point.
As alluded to above, using adaptive routing techniques, a routing methodology is incorporated into a network, such as network 100 (
Again, a switch, such as switch 202, supports minimal and non-minimal path routing based on the destination. If a destination NIC is local, an output port that connects to the destination switch is selected. If the destination is in another group, the packet is routed to a switch within the local group that is connected to the destination group. In a large system, there may only be one such path, but in a small system there are likely to be many, some connected to the input switch and others to switches elsewhere within its group. The input switch selects between them. Candidate ports considered for minimal path routing are further divided into preferred and nonpreferred sets of ports where ports in the preferred set may lead to a path containing fewer hops. Non-minimal paths route packets via an intermediary switch, referred to as a root switch. Root switches are selected on a packet-by-packet basis or a flow-by-flow basis depending on the ordering requirements of the traffic.
A global minimal path between an edge port in one group and an edge port in another group can require one, two, or three hops across fabric links. When identifying minimal path candidate ports, the FRF is able to classify the candidates into a set of preferred ports and a set of non-preferred ports. The preferred ports are those that allow a global minimal path requiring two, or fewer fabric link hops. Use of preferred paths, when available, and when not too heavily loaded reduces the average load on the system's local fabric links as it reduces the average number of local fabric links traversed per frame. Use of a preferred path may also reduce the end-to-end fabric latency experienced by the frame.
It should be noted that the terms “optimize,” “optimal” and the like as used herein can be used to mean making or achieving performance as effective or perfect as possible. However, as one of ordinary skill in the art reading this document will recognize, perfection cannot always be achieved. Accordingly, these terms can also encompass making or achieving performance as good or effective as possible or practical under the given circumstances, or making or achieving performance better than that which can be achieved with other settings or parameters.
The computer system 900 also includes a main memory 906, such as a random access memory (RAM), cache and/or other dynamic storage devices, coupled to bus 902 for storing information and instructions to be executed by processor 904. Main memory 906 also may be used for storing temporary variables or other intermediate information during execution of instructions to be executed by processor 904. Such instructions, when stored in storage media accessible to processor 904, render computer system 900 into a special-purpose machine that is customized to perform the operations specified in the instructions.
The computer system 900 further includes a read only memory (ROM) 908 or other static storage device coupled to bus 902 for storing static information and instructions for processor 904. A storage device 910, such as a magnetic disk, optical disk, or USB thumb drive (Flash drive), etc., is provided and coupled to bus 902 for storing information and instructions.
The computer system 900 may be coupled via bus 902 to a display 912, such as a liquid crystal display (LCD) (or touch screen), for displaying information to a computer user. An input device 914, including alphanumeric and other keys, is coupled to bus 902 for communicating information and command selections to processor 904. Another type of user input device is cursor control 916, such as a mouse, a trackball, or cursor direction keys for communicating direction information and command selections to processor 904 and for controlling cursor movement on display 912. In some embodiments, the same direction information and command selections as cursor control may be implemented via receiving touches on a touch screen without a cursor.
The computing system 900 may include a user interface module to implement a GUI that may be stored in a mass storage device as executable software codes that are executed by the computing device(s). This and other modules may include, by way of example, components, such as software components, object-oriented software components, class components and task components, processes, functions, attributes, procedures, subroutines, segments of program code, drivers, firmware, microcode, circuitry, data, databases, data structures, tables, arrays, and variables.
In general, the word “component,” “engine,” “system,” “database,” data store,” and the like, as used herein, can refer to logic embodied in hardware or firmware, or to a collection of software instructions, possibly having entry and exit points, written in a programming language, such as, for example, Java, C or C++. A software component may be compiled and linked into an executable program, installed in a dynamic link library, or may be written in an interpreted programming language such as, for example, BASIC, Perl, or Python. It will be appreciated that software components may be callable from other components or from themselves, and/or may be invoked in response to detected events or interrupts. Software components configured for execution on computing devices may be provided on a computer readable medium, such as a compact disc, digital video disc, flash drive, magnetic disc, or any other tangible medium, or as a digital download (and may be originally stored in a compressed or installable format that requires installation, decompression or decryption prior to execution). Such software code may be stored, partially or fully, on a memory device of the executing computing device, for execution by the computing device. Software instructions may be embedded in firmware, such as an EPROM. It will be further appreciated that hardware components may be comprised of connected logic units, such as gates and flip-flops, and/or may be comprised of programmable units, such as programmable gate arrays or processors.
The computer system 900 may implement the techniques described herein using customized hard-wired logic, one or more ASICs or FPGAs, firmware and/or program logic which in combination with the computer system causes or programs computer system 900 to be a special-purpose machine. According to one embodiment, the techniques herein are performed by computer system 900 in response to processor(s) 904 executing one or more sequences of one or more instructions contained in main memory 906. Such instructions may be read into main memory 906 from another storage medium, such as storage device 910. Execution of the sequences of instructions contained in main memory 906 causes processor(s) 904 to perform the process steps described herein. In alternative embodiments, hard-wired circuitry may be used in place of or in combination with software instructions.
The term “non-transitory media,” and similar terms, as used herein refers to any media that store data and/or instructions that cause a machine to operate in a specific fashion. Such non-transitory media may comprise non-volatile media and/or volatile media. Non-volatile media includes, for example, optical or magnetic disks, such as storage device 910. Volatile media includes dynamic memory, such as main memory 906. Common forms of non-transitory media include, for example, a floppy disk, a flexible disk, hard disk, solid state drive, magnetic tape, or any other magnetic data storage medium, a CD-ROM, any other optical data storage medium, any physical medium with patterns of holes, a RAM, a PROM, and EPROM, a FLASH-EPROM, NVRAM, any other memory chip or cartridge, and networked versions of the same.
Non-transitory media is distinct from but may be used in conjunction with transmission media. Transmission media participates in transferring information between non-transitory media. For example, transmission media includes coaxial cables, copper wire and fiber optics, including the wires that comprise bus 902. Transmission media can also take the form of acoustic or light waves, such as those generated during radio-wave and infra-red data communications.
The computer system 900 also includes a communication interface 918 coupled to bus 902. Network interface 918 provides a two-way data communication coupling to one or more network links that are connected to one or more local networks. For example, communication interface 918 may be an integrated services digital network (ISDN) card, cable modem, satellite modem, or a modem to provide a data communication connection to a corresponding type of telephone line. As another example, network interface 918 may be a local area network (LAN) card to provide a data communication connection to a compatible LAN (or WAN component to communicated with a WAN). Wireless links may also be implemented. In any such implementation, network interface 918 sends and receives electrical, electromagnetic or optical signals that carry digital data streams representing various types of information.
A network link typically provides data communication through one or more networks to other data devices. For example, a network link may provide a connection through local network to a host computer or to data equipment operated by an Internet Service Provider (ISP). The ISP in turn provides data communication services through the world wide packet data communication network now commonly referred to as the “Internet.” Local network and Internet both use electrical, electromagnetic or optical signals that carry digital data streams. The signals through the various networks and the signals on network link and through communication interface 918, which carry the digital data to and from computer system 900, are example forms of transmission media.
The computer system 900 can send messages and receive data, including program code, through the network(s), network link and communication interface 918. In the Internet example, a server might transmit a requested code for an application program through the Internet, the ISP, the local network and the communication interface 918.
The received code may be executed by processor 904 as it is received, and/or stored in storage device 910, or other non-volatile storage for later execution.
Each of the processes, methods, and algorithms described in the preceding sections may be embodied in, and fully or partially automated by, code components executed by one or more computer systems or computer processors comprising computer hardware. The one or more computer systems or computer processors may also operate to support performance of the relevant operations in a “cloud computing” environment or as a “software as a service” (SaaS). The processes and algorithms may be implemented partially or wholly in application-specific circuitry. The various features and processes described above may be used independently of one another, or may be combined in various ways. Different combinations and sub-combinations are intended to fall within the scope of this disclosure, and certain method or process blocks may be omitted in some implementations. The methods and processes described herein are also not limited to any particular sequence, and the blocks or states relating thereto can be performed in other sequences that are appropriate, or may be performed in parallel, or in some other manner. Blocks or states may be added to or removed from the disclosed example embodiments. The performance of certain of the operations or processes may be distributed among computer systems or computers processors, not only residing within a single machine, but deployed across a number of machines.
As used herein, a circuit might be implemented utilizing any form of hardware, software, or a combination thereof. For example, one or more processors, controllers, ASICs, PLAs, PALs, CPLDs, FPGAs, logical components, software routines or other mechanisms might be implemented to make up a circuit. In implementation, the various circuits described herein might be implemented as discrete circuits or the functions and features described can be shared in part or in total among one or more circuits. Even though various features or elements of functionality may be individually described or claimed as separate circuits, these features and functionality can be shared among one or more common circuits, and such description shall not require or imply that separate circuits are required to implement such features or functionality. Where a circuit is implemented in whole or in part using software, such software can be implemented to operate with a computing or processing system capable of carrying out the functionality described with respect thereto, such as computer system 900.
As used herein, the term “or” may be construed in either an inclusive or exclusive sense. Moreover, the description of resources, operations, or structures in the singular shall not be read to exclude the plural. Conditional language, such as, among others, “can,” “could,” “might,” or “may,” unless specifically stated otherwise, or otherwise understood within the context as used, is generally intended to convey that certain embodiments include, while other embodiments do not include, certain features, elements and/or steps.
Terms and phrases used in this document, and variations thereof, unless otherwise expressly stated, should be construed as open ended as opposed to limiting. Adjectives such as “conventional,” “traditional,” “normal,” “standard,” “known,” and terms of similar meaning should not be construed as limiting the item described to a given time period or to an item available as of a given time, but instead should be read to encompass conventional, traditional, normal, or standard technologies that may be available or known now or at any time in the future. The presence of broadening words and phrases such as “one or more,” “at least,” “but not limited to” or other like phrases in some instances shall not be read to mean that the narrower case is intended or required in instances where such broadening phrases may be absent.
Roweth, Duncan, Froese, Edwin L., Bataineh, Abdulla M.
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